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| CVE | Vendors | Products | Updated | CVSS v3.1 |
|---|---|---|---|---|
| CVE-2026-89808 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: drm/amdkfd: Fix the case that vm range is hole at svm_migrate_copy_to_vram When migration vm range is hole at cpu side(MIGRATE_PFN_MIGRATE set + MIGRATE_PFN_VALID unset) driver still allocates device pages. There is no dma map of src pages and migration. j is 0 and svm_migrate_copy_memory_gart() will return an uninitialized r. That can trigger out_free_vram_pages to drop all VRAM just set up. Initialize r and only call the last svm_migrate_copy_memory_gart if j > 0. Current code postponed the last page to the final copy. This patch flushes on the last page when reach to the end of current drm_buddy_block; avoids another svm_migrate_copy_memory_gart. | ||||
| CVE-2026-89803 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: drm/nouveau: unsubscribe the channel-kill event before the fence context nouveau_channel_del() tears the fence context down first and only drops the channel-kill subscription later, in the middle of the nvif object teardown: if (chan->fence) nouveau_fence(chan->cli->drm)->context_del(chan); ... nvif_object_dtor(&chan->vram); nvif_event_dtor(&chan->kill); The subscribed handler is nouveau_channel_killed(), which calls nouveau_channel_kill() and from there nouveau_fence_context_kill() on chan->fence. A kill event delivered in that window takes fctx->lock and walks fctx->pending on a fence context that context_del() has already freed. Nothing reaches this below Fermi today, because the subscription is gated on FERMI_CHANNEL_GPFIFO and nothing kills a channel there. On Fermi and newer the window is real but narrow, since a kill has to land exactly while the channel is being destroyed. That is reason enough on its own, which is why this carries a Fixes: tag. The last patch in this series subscribes Tesla channels as well; nothing kills those today, so it does not widen the exposure now, but it is the groundwork for a recovery path that would, and the ordering is better fixed before that lands than alongside it. Drop the subscription before anything it depends on is torn down. | ||||
| CVE-2026-89799 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: bpf: Disable preemption in bpf_get_stackid The get_perf_callchain call needs disabled preemption plus we need it disabled as long as we access its returned trace entries buffer. Note the bpf_get_stackid_pe function is executed already with preemption disabled. | ||||
| CVE-2026-89795 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 8.4 High |
| In the Linux kernel, the following vulnerability has been resolved: PCI: Allow per function PCI slots to fix slot reset on s390 On s390 systems, which use a machine level hypervisor, PCI devices are always accessed through a form of PCI pass-through which fundamentally operates on a per PCI function granularity. This is also reflected in the s390 PCI hotplug driver which creates hotplug slots for individual PCI functions. Its reset_slot() function, which is a wrapper for zpci_hot_reset_device(), thus also resets individual functions. Currently, the pci_create_slot() assigns the same pci_slot object to multifunction devices. This approach worked fine on s390 systems that only exposed virtual functions as individual PCI domains to the operating system. Since commit 44510d6fa0c0 ("s390/pci: Handling multifunctions") s390 supports exposing the topology of multifunction PCI devices by grouping them in a shared PCI domain. This creates a problem when resetting a function through the hotplug driver's slot_reset() interface. When attempting to reset a function through the hotplug driver, the shared slot assignment causes the wrong function to be reset instead of the intended one. It also leaks memory as we do create a pci_slot object for the function, but don't correctly free it in pci_slot_release(). Add a flag for struct pci_slot to allow per function PCI slots for functions managed through a hypervisor, which exposes individual PCI functions while retaining the topology. Since we can use all 8 bits for slot 'number' (for ARI devices), change slot 'number' u16 to account for special values PCI_SLOT_PLACEHOLDER and PCI_SLOT_ALL_DEVICES. | ||||
| CVE-2026-89794 | 1 Linux | 1 Linux Kernel | 2026-10-30 | N/A |
| In the Linux kernel, the following vulnerability has been resolved: ksmbd: zero pipe read compound padding Compound response handling extends the last response iov to an eight-byte boundary. smb2_read_pipe() allocates only the payload size, so the alignment padding can expose up to seven bytes of uninitialized kernel heap memory. Allocate the aligned size and clear the unused tail before pinning the response buffer. | ||||
| CVE-2026-89788 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 9.8 Critical |
| In the Linux kernel, the following vulnerability has been resolved: ksmbd: fix tree connection use-after-free in smb2_tree_connect() ksmbd_tree_conn_connect() publishes a new tree connection in sess->tree_conns with a single reference and returns its pointer to smb2_tree_connect(). The handler continues to initialize the object and build the response after publication. A concurrent session logoff can erase the connection and drop that reference, freeing the object while the handler still uses it. BUG: KASAN: slab-use-after-free in smb2_tree_connect+0xe3d/0xf90 smb2_tree_connect (fs/smb/server/smb2pdu.c:2872) handle_ksmbd_work process_one_work worker_thread kthread After xa_store() succeeds, take a second reference before releasing tree_conns_lock. The original reference belongs to the xarray entry and the second belongs to the creating smb2_tree_connect() handler. Keep the references balanced in every path: - On normal exit or an error after publication, smb2_tree_connect() drops its creator reference. Error cleanup also calls ksmbd_tree_conn_disconnect(), which drops the xarray reference only if it removes the exact entry. - SMB2 TREE_DISCONNECT uses the same helper to remove the entry and drop its xarray reference. The request's existing lookup reference remains owned by the request and is released by the existing cleanup. - Session LOGOFF removes each entry and drops its xarray reference. If it wins the race, later cleanup sees that the entry is gone and does not drop that reference again. To enforce this ownership, claim the disconnected state and erase the exact entry atomically under tree_conns_lock. This guarantees one drop for the xarray reference and one drop by each in-flight user, regardless of which teardown path wins. If logoff removes the entry before initialization completes, fail the connect instead of marking the detached object TREE_CONNECTED. | ||||
| CVE-2026-89763 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: KEYS: trusted: Fix TPM teardown ordering trusted_tpm_exit() drops the TPM chip reference and frees the digest array before unregistering the trusted key type. key_type_lookup() holds key_types_sem for reading until the key operation finishes, while unregister_key_type() takes it for writing. It therefore provides the synchronization point that must precede backend teardown. The current order permits this interleaving: CPU 0 CPU 1 trusted_tpm_exit() key_type_lookup("trusted") put_device(&chip->dev) trusted_tpm_seal() kfree(digests) pcrlock() unregister_key_type() tpm_pcr_extend(..., digests) CPU 1 can consequently dereference the freed digest array. The chip can also be released before callbacks stop using it. KASAN reported: BUG: KASAN: slab-use-after-free in tpm_pcr_extend+0x1f0/0x200 Read of size 2 at addr ffff88810872d000 by task poc/89 Call Trace: tpm_pcr_extend+0x1f0/0x200 pcrlock+0x42/0x70 [trusted] trusted_tpm_seal+0x1b6/0x570 [trusted] trusted_instantiate+0x293/0x340 [trusted] __key_instantiate_and_link+0xb2/0x2b0 __key_create_or_update+0x61e/0xb50 __do_sys_add_key+0x1b8/0x310 Allocated by task 88: __kmalloc_noprof+0x1a7/0x490 do_one_initcall+0xa1/0x390 do_init_module+0x2df/0x840 Freed by task 90: kfree+0x131/0x3c0 trusted_tpm_exit+0x59/0xa0 [trusted] __do_sys_delete_module+0x346/0x510 Move unregister_key_type() before releasing either resource. This stops new lookups and waits for in-flight key operations to finish before the backend state is destroyed. | ||||
| CVE-2026-89762 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: apparmor: fix cred UAF caused by begin_current_label_crit_section() AppArmor's begin_current_label_crit_section() is a scary function called from lots of LSM hooks (in particular VFS/socket-related ones) that checks if the label referenced by the current creds is marked FLAG_STALE, and if so, attempts to use aa_replace_current_label() to replace the creds with an updated version that uses a new label. The first problem with this is that it would directly lead to UAF of `struct cred` if anything in the kernel takes a pointer to the current creds and accesses these past a security hook invocation that replaces creds, like so: ``` const struct cred *cred = current_cred(); alloc_file_pseudo(...); uid_t uid = cred->euid; ``` I don't know if anything in the kernel actually does this, but I think it is very surprising that this pattern could lead to UAF. The second problem is that things go wrong when aa_replace_current_label() runs with overridden credentials. aa_replace_current_label() bails out if `current_cred() != current_real_cred()` (mirroring the check in proc_pid_attr_write()), but this check can't actually reliably detect overridden credentials because the overridden creds can be the same as the objective creds. So in approximately the following scenario, things go wrong: 1. task begins with <creds A> (as both objective and subjective creds), with refcount=2 2. task grabs an extra reference on <creds A> for overriding 3. task calls override_creds(<creds A>), which returns a pointer to the old subjective creds (<creds A>) 4. task enters AppArmor LSM hook 5. AppArmor checks that objective/subjective creds are equal 6. AppArmor replaces both cred pointers with <creds B> and drops 2 refs on <creds A> 7. task leaves AppArmor LSM hook 8. task calls revert_creds(<creds A>) 9. now task->cred is <creds A> while task->real_cred is <creds B>, but the task_struct logically holds two references to <creds B> 10. another task drops the extra reference on <creds A> that was used for overriding, refcount drops to 0 11. now task->real_cred points to freed creds At this point, any access to current_cred() will be UAF. I have a test case where I run aa-disable on a profile while a process using that profile is blocked on splice() from a FUSE passthrough file into a full pipe; after the profile update, the pipe becomes empty, splice() resumes, the credentials go out of sync, and a subsequent getuid() syscall results in a KASAN UAF splat. To fix this, instead of directly replacing creds, do it via task_work that will run at the end of the current syscall. (The point in time at which the cred replacement happens should have no correctness impact; it is just a performance optimization to avoid unnecessarily touching the refcount of the new label.) Note that AppArmor still performs direct cred replacements in the sb_pivotroot LSM hook after this change, and that direct cred replacements can still happen in VFS ->write() callbacks via proc_pid_attr_write(). There are two options for what to do with aa_dup_task_ctx(): Either explicitly reset new->label_replacement_pending after the entire aa_task_ctx has been copied, or switch to manually copying members over. I am switching to manually copying members over because that should make bugs more obvious. | ||||
| CVE-2026-89761 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: apparmor: fix out-of-bounds write when null terminating a label vec aa_vec_unique() null terminates at vec[n - dups] when VEC_FLAG_TERMINATE is passed. If the components are all distinct no duplicates are dropped, dups is 0 and the terminator goes to vec[n], so the caller has to provide room for n + 1 entries. aa_label_strn_parse() sets up its vector with vec_setup(profile, vec, len, gfp) and then calls aa_vec_unique(vec, len, VEC_FLAG_TERMINATE), but vec_setup() does not reserve the terminator entry. Up to LOCAL_VEC_ENTRIES it uses the local array of LOCAL_VEC_ENTRIES pointers, above that it allocates exactly len pointers. The terminator therefore lands one entry past the end of the local array when len is LOCAL_VEC_ENTRIES, and one entry past the end of the allocation when len is larger. len comes from the number of "//&" separated components in the label name and label_count_strn_entries() does not bound it. An unprivileged task reaches the parse by writing to /proc/self/attr/apparmor/current or through lsm_set_self_attr(2), both of which go through do_setattr(), and the name is parsed before the change_profile permission is checked. The query_label() path behind the securityfs .access file, which is mode 0666, performs no permission check at all. Every component has to resolve to a loaded profile, so a system with policy loaded is required. The other two VEC_FLAG_TERMINATE users work on a label vec that aa_label_alloc() has already sized with "+ 1 for null terminator entry on vec". Reserve the same entry in vec_setup() and DEFINE_VEC(). Passing len + 1 from the caller instead would move len == LOCAL_VEC_ENTRIES out of the local array and into kzalloc(). | ||||
| CVE-2026-89759 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 4.7 Medium |
| In the Linux kernel, the following vulnerability has been resolved: mm/kmemleak: avoid soft lockup when scanning task stacks Patch series "mm/kmemleak: avoid soft lockup when scanning task", v3. kmemleak_scan() scans every task stack under one rcu_read_lock() with no reschedule point, which can trip the soft lockup watchdog on hosts with very many threads. That prints the following message, depending on the workload+host configuration: watchdog: BUG: soft lockup - CPU#35 stuck for 22s! [kmemleak:537] scan_block kmemleak_scan kmemleak_scan_thread kthread Patch 1 walks the tasks with find_ge_pid() so the scan reschedules between tasks Patches 2-3 let the scan loops stop early once a scan is interrupted. This patch (of 3): kmemleak_scan() walks every thread and scans its kernel stack under a single rcu_read_lock() with no reschedule point. On a host with very many threads -- amplified by KASAN/lockdep in debug builds -- this loop can hog a CPU long enough to trip the soft lockup watchdog: watchdog: BUG: soft lockup - CPU#35 stuck for 22s! [kmemleak:537] scan_block kmemleak_scan kmemleak_scan_thread kthread A cond_resched() cannot be added directly: the loop runs inside an RCU read-side critical section. Walk the tasks one PID at a time with find_ge_pid(), taking the RCU read lock only to look up and pin each task. The stack is then scanned with no lock held, so cond_resched() runs between tasks and the scan stops early on scan_should_stop(). This follows the next_tgid()/task_seq_get_next() iteration pattern and keeps each RCU critical section short. | ||||
| CVE-2026-89757 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 5.3 Medium |
| In the Linux kernel, the following vulnerability has been resolved: mm/mglru: fix and remove redundant unevictable folio handling sort_folio() has a shortcut for moving folios that are no longer evictable but are still sitting on a generation list. However, this shortcut is buggy. It does not follow the PG_lru usage convention, and it has a more serious issue. Unevictable folios are not threaded on lists[LRU_UNEVICTABLE], so that folio->lru can be reused to hold folio->mlock_count (see the comment in lruvec_init()). Hence lruvec_add_folio() skips the list_add() for them, and every other place that turns a folio unevictable initialises mlock_count explicitly: lru_add() sets it to 0, __mlock_folio() and __mlock_new_folio() set it to !!folio_test_mlocked(folio). sort_folio() sets nothing, and the lru_gen_del_folio() right above it may have already poisoned folio->lru via list_del(), so mlock_count ends up aliasing LIST_POISON2, which reads as 0x122, i.e. 290. The result is user visible. On munlock, __munlock_folio() decrements that bogus count, finds it still non-zero and bails out before clearing PG_mlocked, so the folio remains unevictable and the Mlocked accounting stays inflated until the folio is freed. The shortcut also touches the LRU flags in the wrong order. It calls lru_gen_del_folio() while PG_lru is still set, so a concurrent folio_test_clear_lru() (e.g. compaction, folio_isolate_lru()) can succeed on a folio that has already been taken off the generation list, which may lead to unexpected behavior. So fix it by isolating them as common folios and letting the generic shrink path cull them. This matches the classical LRU behavior, and there should be no visible effect on the generic eviction or isolation behavior. There is no performance concern either, such a folio goes through this once, and then it is off the generation lists for good. | ||||
| CVE-2026-89744 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 8.4 High |
| In the Linux kernel, the following vulnerability has been resolved: device property: fix infinite loop in fwnode_for_each_child_node() When iterate over children of a fwnode that has a secondary fwnode, fwnode_get_next_child_node() can enter an infinite loop if the secondary fwnode has more than one child. Parent Child (Primary fwnode) FWa: {FWa1, FWa2, FWa3} (Secondary fwnode) FWb: {FWb1, FWb2} In this case: ┌─> fwnode_get_next_child_node(FWa, FWa1) │ - fwnode_call_ptr_op(FWa, get_next_child_node, FWa1) returns FWa2 │ │ ... │ │ fwnode_get_next_child_node(FWa, FWa3) │ - fwnode_call_ptr_op(FWa, get_next_child_node, FWa3) returns NULL │ - fwnode_call_ptr_op(FWb, get_next_child_node, FWa3) returns FWb1 │ │ fwnode_get_next_child_node(FWa, FWb1) │ - fwnode_call_ptr_op(FWa, get_next_child_node, FWb1) returns FWa1 └────┘ This cause fwnode_for_each_child_node() to loop indefinitely, reapeatedly output {FWa1, FWa2, FWa3, FWb1, FWa1, ...}. The root cause is that when the current child (FWb1) belongs to the secondary fwnode, calling get_next_child_node() on the parimary fwnode incorrectly returns the first child (FWa1) again instead of NULL. Fix this by dynamically checking the parent fwnode of the current child before calling get_next_child_node(). This approach follows the pattern established in commit b5b41ab6b0c1 ("device property: Check fwnode->secondary in fwnode_graph_get_next_endpoint()"). | ||||
| CVE-2026-89740 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 6.3 Medium |
| In the Linux kernel, the following vulnerability has been resolved: serial: imx: serialize imx_uart_ports[] lifetime imx_uart_probe() publishes its devm-allocated port in imx_uart_ports[] before uart_add_one_port() because console setup uses the table. The entry is not cleared when adding the port fails or after removal, leaving a dangling pointer. A sibling probe can register the shared console through that stale entry. This was reproduced under KASAN on QEMU mcimx6ul-evk by unbinding a sibling UART, unbinding the console UART and rebinding the sibling. Keep the entry valid through uart_remove_one_port(), then clear it. Protect port addition and removal together with their table updates so sibling operations cannot interleave. Reject an occupied slot rather than clobbering an active port during a duplicate-line probe. | ||||
| CVE-2026-89738 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 7.8 High |
| In the Linux kernel, the following vulnerability has been resolved: usb: gadget: at91_udc: drain polled-VBUS timer/work before udc is freed In polled-VBUS mode (board.vbus_pin && board.vbus_polled), probe arms a self-restarting cycle: at91_vbus_timer() schedules vbus_timer_work, and at91_vbus_timer_work() calls at91_vbus_update() and re-arms the timer via mod_timer(). Both recover the same udc through container_of and dereference it on every iteration. Neither teardown path cancels this cycle. udc is devm-allocated, so it is freed after at91udc_remove() returns, and is likewise freed when probe fails and devres runs. A timer callback or work item that is pending or running at either point dereferences the freed udc. Add at91_udc_shutdown_vbus_timer() and call it from at91udc_remove() and from the usb_add_gadget_udc() failure path in probe; the remaining probe error paths fail before the timer is armed. timer_shutdown_sync() waits for a running callback and clears timer->function, which makes the work handler's mod_timer() a permanent no-op; cancel_work_sync() then drains any pending or running work whose re-arm attempt now does nothing. The timer must be shut down first, since cancelling the work alone would let the timer re-queue it. The guard mirrors probe: in IRQ mode the timer and work_struct are never initialized. This does not require a fault; a normal driver unbind can interleave with an already queued work item. This issue was found by an in-house static analysis tool. | ||||
| CVE-2026-89719 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 4.1 Medium |
| In the Linux kernel, the following vulnerability has been resolved: zram: fix out-of-bounds access in read_block_state() read_block_state() calculates nr_pages before taking dev_lock. If the device is reset and reinitialized with a smaller disksize before lock acquisition, nr_pages still describes the old table. The subsequent loop can then call slot_lock() past the end of the newly allocated table. Read disksize after acquiring dev_lock and checking that the device is initialized. The read lock then keeps the table and its bound stable for the duration of the scan. | ||||
| CVE-2026-89718 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 4.1 Medium |
| In the Linux kernel, the following vulnerability has been resolved: zram: fix out-of-bounds access in writeback_store() Patch series "zram: fix stale scan bounds after reinitialization". Both writeback_store() and read_block_state() derive their table scan bounds from zram->disksize before acquiring dev_lock. If the device is reset and reinitialized with a smaller disksize between that read and lock acquisition, the bound can describe the old table while the scan operates on the new one. This can lead to out-of-bounds slot accesses. Move both bound calculations under dev_lock so each bound remains consistent with the table throughout its scan. Keep the fixes separate because the affected interfaces originate from different commits and can be backported independently. This patch (of 2): writeback_store() calculates the table scan bounds before taking dev_lock. A reset followed by reconfiguration with a smaller disksize can therefore replace zram->table while writeback_store() is waiting for the lock. Once it acquires the lock, it sees an initialized device but scans the new table using the old upper bound, resulting in an out-of-bounds access. Calculate the number of pages while holding dev_lock so the scan bound matches the table protected by the lock. | ||||
| CVE-2026-89713 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 9.1 Critical |
| In the Linux kernel, the following vulnerability has been resolved: NFSD: check truncate permission under inode lock nfsd_setattr() checks whether a size update needs NFSD_MAY_TRUNC before it takes inode_lock(). The comparison uses the file size sampled by that unlocked read, but the actual ATTR_SIZE update is applied later under inode_lock() by notify_change(). This leaves a TOCTOU window for append-only files. If a client sends a SETATTR that does not shrink the file at the time of the unlocked sample, a concurrent append can extend the file before nfsd_setattr() takes inode_lock(). notify_change() then applies a real truncation without the NFSD_MAY_TRUNC check that rejects IS_APPEND(inode). The VFS truncate syscall paths perform their own append-only checks before calling notify_change(), so NFSD must make this decision against the locked size it is about to change. Split the write-count acquisition from the truncation permission check. Keep get_write_access() before the locked setattr work, then recheck whether the requested size is below i_size_read(inode) after inode_lock() has been acquired and before notify_change(ATTR_SIZE). This also avoids the plain unlocked inode->i_size load. | ||||
| CVE-2026-89709 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 8.1 High |
| In the Linux kernel, the following vulnerability has been resolved: lockd, nfsd: RCU-protect nlmsvc_ops dispatch nlmsvc_ops is published by nfsd_lockd_init() and cleared by nfsd_lockd_shutdown() with plain stores, while lockd dereferences it unguarded from dispatch sites in fs/lockd/svcsubs.c. The pointer targets nfsd's .rodata and the fopen/fclose callbacks live in nfsd's .text, so a stale load after rmmod nfsd results in either a NULL deref or a module-text use-after-free. Declare nlmsvc_ops as __rcu, publish via rcu_assign_pointer(), clear via RCU_INIT_POINTER() + synchronize_rcu(). Add a struct module *owner field to nlmsvc_binding and pin the module across indirect calls with try_module_get/module_put. When the binding is torn down, fall back to fput() to avoid leaking struct file references. | ||||
| CVE-2026-89707 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 7.5 High |
| In the Linux kernel, the following vulnerability has been resolved: nfsd: release path refs on follow_down() error nfsd_cross_mnt() initializes a local struct path with mntget() and dget() before calling follow_down(). On a negative return the error arm jumps to out without releasing those references: err = follow_down(&path, follow_flags); if (err < 0) goto out; follow_down() never drops the caller's entry-time refs on any error sub-case; for example a pre-cross d_manage() failure leaves path untouched, so the mntget()/dget() taken on entry survive the call. Every other early-exit arm in nfsd_cross_mnt() (other-namespace return, IS_ERR(exp2), and the success tail after the swap) already calls path_put(&path); the err < 0 arm is the lone omission. The leak inflates mnt_count and d_count on each failed cross-mount, blocking umount and pinning dentries against the shrinker, and is reachable by any authenticated NFS client through nfsd_lookup_dentry or the NFSv4 READDIR encode path. Fix by calling path_put(&path) before the goto out in the err < 0 arm so the entry-time refs are released on all follow_down() error returns. | ||||
| CVE-2026-89702 | 1 Linux | 1 Linux Kernel | 2026-10-30 | 9.8 Critical |
| In the Linux kernel, the following vulnerability has been resolved: nfsd: size fh_verify server sockaddr slot by xpt_locallen The nfsd_fh_verify and nfsd_fh_verify_err tracepoints declare the server sockaddr slot sized by xpt_remotelen but fill it from xpt_local using xpt_locallen: TP_STRUCT__entry( ... __sockaddr(server, rqstp->rq_xprt->xpt_remotelen) ... ) TP_fast_assign( ... __assign_sockaddr(server, &rqstp->rq_xprt->xpt_local, rqstp->rq_xprt->xpt_locallen); ... ) When xpt_locallen exceeds xpt_remotelen, __assign_sockaddr's memcpy writes past the reserved ring-buffer slot. In the reverse direction (xpt_locallen < xpt_remotelen) the slot is oversized and the unwritten tail leaks prior ring-buffer contents to trace consumers. The write-past-end case is reachable on NFS/UDP. svc_xprt_set_remote() is only called from svc_tcp_accept() (net/sunrpc/svcsock.c) and from the RDMA connect path; svc_create_socket() for UDP calls only svc_xprt_set_local(), so xpt_remotelen stays 0 for the xprt's lifetime. Every fh_verify trace for an NFSv2/v3-over-UDP request then copies 16 or 28 bytes from xpt_local into a zero-byte slot. The other NFSD tracepoints that record the server address (NFSD_TRACE_PROC_CALL_FIELDS, NFSD_TRACE_PROC_RES_FIELDS, SVC_RQST_ENDPOINT_FIELDS) already size the server slot by xpt_locallen; nfsd_fh_verify and nfsd_fh_verify_err were the only exceptions. Fix by sizing the server slot with xpt_locallen so the declared slot matches the copy length. The client slot and its assignment already agree on xpt_remotelen and are left untouched. | ||||